Technique for preserving efficiency for replication between clusters of a network

ABSTRACT

A technique preserves efficiency for replication of data between a source node of a source cluster (“source”) and a destination node of a destination cluster (“destination”) of a clustered network. Replication in the clustered network may be effected by leveraging global in-line deduplication at the source to identify and avoid copying duplicate data from the source to the destination. To ensure that the copy of the data on the destination is synchronized with the data received at the source, the source creates a snapshot of the data for use as a baseline copy at the destination. Thereafter, new data received at the source that differs from the baseline snapshot are transmitted and copied to the destination. In addition, the source and destination nodes negotiate to establish a mapping of name-to-data when transferring data (i.e., an extent) between the clusters. Illustratively, the name is an extent key for the extent, such that the negotiated mapping established by the source and destination is based on the extent key associated with the extent.

RELATED APPLICATION

The present application claims priority from commonly owned ProvisionalPatent Application No. 62/199,408, entitled TECHNIQUE FOR PRESERVINGEFFICIENCY FOR REPLICATION BETWEEN CLUSTERS OF A NETWORK, filed on Jul.31, 2015, the contents of which are incorporated herein by reference.

BACKGROUND

Technical Field

The present disclosure relates to replication of data between storagesystems and, more specifically, to data replication between storagesystems of different clusters.

Background Information

A plurality of storage systems may be interconnected as a cluster andconfigured to provide storage service relating to the organization ofdata stored on storage devices coupled to the systems. Each storagesystem may implement a high-level module, such as a file system, tologically organize the data stored on the devices as storage containers,such as volumes and/or logical units (LUNs). The storage systems of thecluster may cooperate to further provide a global-deduplication filesystem.

To improve reliability and facilitate recovery in the event of a failureof the storage system, some or all of the underlying data of the filesystem may be replicated (copied) to another storage system. Forexample, a source storage system may create a restorable image of itsfile system and transmit a copy of that image over a network to adestination storage system. The image may be updated with changesreplicated to the destination storage system. However, duplicate data inthe update may still be replicated to the destination storage system(i.e., data already at the destination may be included in the update).It is desirable to improve replication efficiency by not sendingduplicate data when updating an image previously transmitted betweensource and destination storage systems.

BRIEF DESCRIPTION OF THE DRAWINGS

The above and further advantages of the embodiments herein may be betterunderstood by referring to the following description in conjunction withthe accompanying drawings in which like reference numerals indicateidentically or functionally similar elements, of which:

FIG. 1 is a block diagram of a plurality of nodes interconnected as acluster;

FIG. 2 is a block diagram of a node;

FIG. 3 is a block diagram of a storage input/output (I/O) stack of thenode;

FIG. 4 illustrates a write path of the storage I/O stack;

FIG. 5 illustrates a read path of the storage I/O stack;

FIG. 6 is a block diagram of a volume metadata entry;

FIG. 7 is a block diagram of a dense tree metadata structure;

FIG. 8 is a block diagram of a top level of the dense tree metadatastructure;

FIG. 9 illustrates mapping between levels of the dense tree metadatastructure;

FIG. 10 illustrates a workflow for inserting a volume metadata entryinto the dense tree metadata structure in accordance with a writerequest;

FIG. 11 illustrates merging between levels of the dense tree metadatastructure;

FIG. 12 is a block diagram of a dense tree metadata structure sharedbetween a parent volume and snapshot/clone;

FIG. 13 illustrates diverging of the snapshot/clone from the parentvolume; and

FIG. 14 is a block diagram of a technique for preserving efficiency ofreplication between a source cluster and destination cluster of aclustered network.

OVERVIEW

The embodiments herein are directed to a technique for preservingefficiency for replication of data between a source node of a sourcecluster (“source”) and a destination node of a destination cluster(“destination”) of a clustered network. Data replication in theclustered network may be performed by leveraging global deduplication ofthe cluster to identify and avoid copying duplicate data from the sourceto the destination. To ensure that the copy of the data on thedestination is synchronized with the data received at the source, thesource may create a snapshot of the data for use as a baseline copy atthe destination. Thereafter, new data received at the source thatdiffers from the baseline snapshot (i.e., copy) are transmitted andcopied to the destination. The new data may be data associated with oneor more write requests (i.e., write data) issued by a host and directedto a logical block address range of a logical unit served by the source.The write data may be organized, e.g. aggregated, into one or morevariable length blocks or extents, which may be de-duplicated.Illustratively, a hash function may be applied to each extent togenerate an extent key that is stored in a key-value extent store (ES)embodied as a data structure, e.g., an ES hash table, of each cluster.The extent key is configured to reference a location of the extent onone or more storage devices, such as solid state drives of the cluster.As such, replication may occur between two different extent stores ondifferent (e.g., source and destination) clusters, each using the sameextent keys, i.e., a same hash function is used on both clusters.

To preserve efficiency during data replication, the nodes of theclusters may negotiate (e.g., during an initialization stage ofreplication) to ensure that the same hash function is used by the sourceand destination. In addition, the source and destination nodes maynegotiate to establish a mapping of name-to-data when transferring data(i.e., an extent) between the clusters. Illustratively, the name is theextent key for the extent, such that the negotiated mapping establishedby the source and destination is based on the extent key associated withthe extent. To avoid name collisions, the source sends the extent alongwith the extent key (i.e., name) to the destination for the firsttransfer of new data to verify the association (i.e., mapping) of thekey to the extent. The destination accepts the mapping if it can use theextent key-to-extent association (i.e., as a new mapping or as aduplicate of an existing mapping). The mapping may be considered validand in effect when the source and destination agree on the association,and may be considered invalid when, e.g., the extent is deleted byeither the source or destination. The extent is considered duplicatewhen there is an existing mapping associated with the extent key of theextent.

In an embodiment, a replication field is provided within each entry ofthe ES hash table, wherein the replication field is associated on aper-cluster pair basis, e.g., between the source and destinationclusters. The replication field of the ES hash table may include one ormore replication bits organized as a bit plane, e.g., one byte (8 bits)per entry of the ES hash table, wherein each bit represents aper-cluster pair replication relationship, e.g., up to 8 replicationrelationships may be represented using a one byte replication field.Illustratively, each replication bit of the bit plane on the source islinked to a particular destination cluster, which may be indicated usingan associated cluster identifier (ID), e.g., a first replication bit maycorrespond to destination cluster ID X and a second replication bit maycorrespond to destination cluster ID Y. That is, the first replicationbit represents a “Source-X” per-cluster pair replication relationship,whereas the second replication bit represents a “Source-Y” per-clusterpair replication relationship. Accordingly, when the correspondingreplication bits are asserted (e.g., set) for a given extent key in eachES hash table (i.e., in the source ES hash table and in the destinationES hash table), the source and destination agree that the same extentkey is used for the same extent between the clusters. Any otherarrangement of the replication bits (e.g., at least one bit unasserted)requires renegotiation between the source and the destination toestablish the mapping of an extent key to the extent.

DESCRIPTION Storage Cluster

FIG. 1 is a block diagram of a plurality of nodes 200 interconnected asa cluster 100 and configured to provide storage service relating to theorganization of information on storage devices. The nodes 200 may beinterconnected by a cluster interconnect fabric 110 and includefunctional components that cooperate to provide a distributed storagearchitecture of the cluster 100, which may be deployed in a storage areanetwork (SAN). As described herein, the components of each node 200include hardware and software functionality that enable the node toconnect to one or more hosts 120 over a computer network 130, as well asto one or more storage arrays 150 of storage devices over a storageinterconnect 140, to thereby render the storage service in accordancewith the distributed storage architecture.

Each host 120 may be embodied as a general-purpose computer configuredto interact with any node 200 in accordance with a client/server modelof information delivery. That is, the client (host) may request theservices of the node, and the node may return the results of theservices requested by the host, by exchanging packets over the network130. The host may issue packets including file-based access protocols,such as the Network File System (NFS) protocol over the TransmissionControl Protocol/Internet Protocol (TCP/IP), when accessing informationon the node in the form of storage containers such as files anddirectories. However, in an embodiment, the host 120 illustrativelyissues packets including block-based access protocols, such as the SmallComputer Systems Interface (SCSI) protocol encapsulated over TCP (iSCSI)and SCSI encapsulated over FC (FCP), when accessing information in theform of storage containers such as logical units (LUNs). Notably, any ofthe nodes 200 may service a request directed to a storage containerstored on the cluster 100.

FIG. 2 is a block diagram of a node 200 that is illustratively embodiedas a storage system having one or more central processing units (CPUs)210 coupled to a memory 220 via a memory bus 215. The CPU 210 is alsocoupled to a network adapter 230, storage controllers 240, a clusterinterconnect interface 250 and a non-volatile random access memory(NVRAM 280) via a system interconnect 270. The network adapter 230 mayinclude one or more ports adapted to couple the node 200 to the host(s)120 over computer network 130, which may include point-to-point links,wide area networks, virtual private networks implemented over a publicnetwork (Internet) or a local area network. The network adapter 230 thusincludes the mechanical, electrical and signaling circuitry needed toconnect the node to the network 130, which illustratively embodies anEthernet or Fibre Channel (FC) network.

The memory 220 may include memory locations that are addressable by theCPU 210 for storing software programs and data structures associatedwith the embodiments described herein. The CPU 210 may, in turn, includeprocessing elements and/or logic circuitry configured to execute thesoftware programs, such as a storage input/output (I/O) stack 300, andmanipulate the data structures. Illustratively, the storage I/O stack300 may be implemented as a set of user mode processes that may bedecomposed into a plurality of threads. An operating system kernel 224,portions of which are typically resident in memory 220 (in-core) andexecuted by the processing elements (i.e., CPU 210), functionallyorganizes the node by, inter alia, invoking operations in support of thestorage service implemented by the node and, in particular, the storageI/O stack 300. A suitable operating system kernel 224 may include ageneral-purpose operating system, such as the UNIX® series or MicrosoftWindows® series of operating systems, or an operating system withconfigurable functionality such as microkernels and embedded kernels.However, in an embodiment described herein, the operating system kernelis illustratively the Linux® operating system. It will be apparent tothose skilled in the art that other processing and memory means,including various computer readable media, may be used to store andexecute program instructions pertaining to the embodiments herein.

Each storage controller 240 cooperates with the storage I/O stack 300executing on the node 200 to access information requested by the host120. The information is preferably stored on storage devices such assolid state drives (SSDs) 260, illustratively embodied as flash storagedevices, of storage array 150. In an embodiment, the flash storagedevices may be based on NAND flash components, e.g., single-layer-cell(SLC) flash, multi-layer-cell (MLC) flash or triple-layer-cell (TLC)flash, although it will be understood to those skilled in the art thatother non-volatile, solid-state electronic devices (e.g., drives basedon storage class memory components) may be advantageously used with theembodiments described herein. Accordingly, the storage devices may ormay not be block-oriented (i.e., accessed as blocks). The storagecontroller 240 includes one or more ports having I/O interface circuitrythat couples to the SSDs 260 over the storage interconnect 140,illustratively embodied as a serial attached SCSI (SAS) topology.Alternatively, other point-to-point I/O interconnect arrangements, suchas a conventional serial ATA (SATA) topology or a PCI topology, may beused. The system interconnect 270 may also couple the node 200 to alocal service storage device 248, such as an SSD, configured to locallystore cluster-related configuration information, e.g., as clusterdatabase (DB) 244, which may be replicated to the other nodes 200 in thecluster 100.

The cluster interconnect interface 250 may include one or more portsadapted to couple the node 200 to the other node(s) of the cluster 100.In an embodiment, Infiniband may be used as the clustering protocol andinterconnect fabric media, although it will be apparent to those skilledin the art that other types of protocols and interconnects may beutilized within the embodiments described herein. The NVRAM 280 mayinclude a back-up battery or other built-in last-state retentioncapability (e.g., non-volatile semiconductor memory such as storageclass memory) that is capable of maintaining data in light of a failureto the node and cluster environment. Illustratively, a portion of theNVRAM 280 may be configured as one or more non-volatile logs (NVLogs285) configured to temporarily record (“log”) I/O requests, such aswrite requests, received from the host 120.

Storage I/O Stack

FIG. 3 is a block diagram of the storage I/O stack 300 that may beadvantageously used with one or more embodiments described herein. Thestorage I/O stack 300 includes a plurality of software modules or layersthat cooperate with other functional components of the nodes 200 toprovide the distributed storage architecture of the cluster 100. In anembodiment, the distributed storage architecture presents an abstractionof a single storage container, i.e., all of the storage arrays 150 ofthe nodes 200 for the entire cluster 100 organized as one large pool ofstorage. In other words, the architecture consolidates storage, i.e.,the SSDs 260 of the arrays 150, throughout the cluster (retrievable viacluster-wide keys) to enable storage of the LUNs. Both storage capacityand performance may then be subsequently scaled by adding nodes 200 tothe cluster 100.

Illustratively, the storage I/O stack 300 includes an administrationlayer 310, a replication layer 315 (described later herein), a protocollayer 320, a persistence layer 330, a volume layer 340, an extent storelayer 350, a Redundant Array of Independent Disks (RAID) layer 360, astorage layer 365 and a NVRAM (storing NVLogs) “layer” interconnectedwith a messaging kernel 370. The messaging kernel 370 may provide amessage-based (or event-based) scheduling model (e.g., asynchronousscheduling) that employs messages as fundamental units of work exchanged(i.e., passed) among the layers. Suitable message-passing mechanismsprovided by the messaging kernel to transfer information between thelayers of the storage I/O stack 300 may include, e.g., for intra-nodecommunication: i) messages that execute on a pool of threads, ii)messages that execute on a single thread progressing as an operationthrough the storage I/O stack, iii) messages using an Inter ProcessCommunication (IPC) mechanism and, e.g., for inter-node communication:messages using a Remote Procedure Call (RPC) mechanism in accordancewith a function shipping implementation. Alternatively, the I/O stackmay be implemented using a thread-based or stack-based execution model.In one or more embodiments, the messaging kernel 370 allocatesprocessing resources from the operating system kernel 224 to execute themessages. Each storage I/O stack layer may be implemented as one or moreinstances (i.e., processes) executing one or more threads (e.g., inkernel or user space) that process the messages passed between thelayers such that the messages provide synchronization for blocking andnon-blocking operation of the layers. Note that one or more of thelayers, such as the administrative layer 310 and the replication layer315, may execute asynchronously to the other layers.

In an embodiment, the protocol layer 320 may communicate with the host120 over the network 130 by exchanging discrete frames or packetsconfigured as I/O requests according to pre-defined protocols, such asiSCSI and FCP. An I/O request, e.g., a read or write request, may bedirected to a LUN and may include I/O parameters such as, inter alia, aLUN identifier (ID), a logical block address (LBA) of the LUN, a length(i.e., amount of data) and, in the case of a write request, write data.The protocol layer 320 receives the I/O request and forwards it to thepersistence layer 330, which records the request into a persistentwrite-back cache 380, illustratively embodied as a log whose contentscan be replaced randomly, e.g., under some random access replacementpolicy rather than only in log fashion, and returns an acknowledgementto the host 120 via the protocol layer 320. In an embodiment only I/Orequests that modify the LUN, e.g., write requests, are logged. Notably,the I/O request may be logged at the node receiving the I/O request, orin an alternative embodiment in accordance with the function shippingimplementation, the I/O request may be logged at another node.

Illustratively, dedicated logs may be maintained by the various layersof the storage I/O stack 300. For example, a dedicated log 335 may bemaintained by the persistence layer 330 to record the I/O parameters ofan I/O request as equivalent internal, i.e., storage I/O stack,parameters, e.g., volume ID, offset, and length. In the case of a writerequest, the persistence layer 330 may also cooperate with the NVRAM 280to implement the write-back cache 380 configured to store the write dataassociated with the write request. Notably, the write data for the writerequest may be physically stored in the log 355 such that the cache 380contains the reference to the associated write data. That is, thewrite-back cache may be structured as a log. In an embodiment, a copy ofthe write-back cache may be also maintained in the memory 220 tofacilitate direct memory access to the storage controllers. In otherembodiments, caching may be performed at the host 120 or at a receivingnode in accordance with a protocol that maintains coherency between thewrite data stored at the cache and the cluster.

In an embodiment, the administration layer 310 may apportion the LUNinto multiple volumes, each of which may be partitioned into multipleregions (e.g., allotted as disjoint block address ranges), with eachregion having one or more segments stored as multiple stripes on thearray 150. A plurality of volumes distributed among the nodes 200 maythus service a single LUN, i.e., each volume within the LUN services adifferent LBA range (i.e., offset and length, hereinafter offset andrange) or set of ranges within the LUN. Accordingly, the protocol layer320 may implement a volume mapping technique to identify a volume towhich the I/O request is directed (i.e., the volume servicing the offsetrange indicated by the parameters of the I/O request). Illustratively,the cluster database 244 may be configured to maintain one or moreassociations (e.g., key-value pairs) for each of the multiple volumes,e.g., an association between the LUN ID and a volume, as well as anassociation between the volume and a node ID for a node managing thevolume. The administration layer 310 may also cooperate with thedatabase 244 to create (or delete) one or more volumes associated withthe LUN (e.g., creating a volume ID/LUN key-value pair in the database244). Using the LUN ID and LBA (or LBA range), the volume mappingtechnique may provide a volume ID (e.g., using appropriate associationsin the cluster database 244) that identifies the volume and nodeservicing the volume destined for the request, as well as translate theLBA (or LBA range) into an offset and length within the volume.Specifically, the volume ID is used to determine a volume layer instancethat manages volume metadata associated with the LBA or LBA range. Asnoted, the protocol layer may pass the I/O request (i.e., volume ID,offset and length) to the persistence layer 330, which may use thefunction shipping (e.g., inter-node) implementation to forward the I/Orequest to the appropriate volume layer instance executing on a node inthe cluster based on the volume ID.

In an embodiment, the volume layer 340 may manage the volume metadataby, e.g., maintaining states of host-visible containers, such as rangesof LUNs, and performing data management functions, such as creation ofsnapshots and clones, for the LUNs in cooperation with theadministration layer 310. The volume metadata is illustratively embodiedas in-core mappings from LUN addresses (i.e., LBAs) to durable extentkeys, which are unique cluster-wide IDs associated with SSD storagelocations for extents within an extent key space of the cluster-widestorage container. That is, an extent key may be used to retrieve thedata of the extent at an SSD storage location associated with the extentkey. Alternatively, there may be multiple storage containers in thecluster wherein each container has its own extent key space, e.g., wherethe host provides distribution of extents among the storage containersand cluster-wide (across containers) de-duplication is infrequent. Anextent is a variable length block of data that provides a unit ofstorage on the SSDs and that need not be aligned on any specificboundary, i.e., it may be byte aligned. Accordingly, an extent may be anaggregation of write data from a plurality of write requests to maintainsuch alignment. Illustratively, the volume layer 340 may record theforwarded request (e.g., information or parameters characterizing therequest), as well as changes to the volume metadata, in dedicated log345 maintained by the volume layer 340. Subsequently, the contents ofthe volume layer log 345 may be written to the storage array 150 inaccordance with retirement of log entries, while a checkpoint (e.g.,synchronization) operation stores in-core metadata on the array 150.That is, the checkpoint operation (checkpoint) ensures that a consistentstate of metadata, as processed in-core, is committed to (stored on) thestorage array 150; whereas the retirement of log entries ensures thatthe entries accumulated in the volume layer log 345 synchronize with themetadata checkpoints committed to the storage array 150 by, e.g.,retiring those accumulated log entries prior to the checkpoint. In oneor more embodiments, the checkpoint and retirement of log entries may bedata driven, periodic or both.

In an embodiment, the extent store layer 350 is responsible for storingextents on the SSDs 260 (i.e., on the storage array 150) and forproviding the extent keys to the volume layer 340 (e.g., in response toa forwarded write request). The extent store layer 350 is alsoresponsible for retrieving data (e.g., an existing extent) using anextent key (e.g., in response to a forwarded read request). In analternative embodiment, the extent store layer 350 is responsible forperforming de-duplication and compression on the extents prior tostorage. The extent store layer 350 may maintain in-core mappings (e.g.,embodied as hash tables) of extent keys to SSD storage locations (e.g.,offset on an SSD 260 of array 150). The extent store layer 350 may alsomaintain a dedicated log 355 of entries that accumulate requested “put”and “delete” operations (i.e., write requests and delete requests forextents issued from other layers to the extent store layer 350), wherethese operations change the in-core mappings (i.e., hash table entries).Subsequently, the in-core mappings and contents of the extent storelayer log 355 may be written to the storage array 150 in accordance witha “fuzzy” checkpoint 390 (i.e., checkpoint with incremental changes thatspan multiple log files) in which selected in-core mappings, less thanthe total, are committed to the array 150 at various intervals (e.g.,driven by an amount of change to the in-core mappings, size thresholdsof log 355, or periodically). Notably, the accumulated entries in log355 may be retired once all in-core mappings have been committed andthen, illustratively, for those entries prior to the first interval.

In an embodiment, the RAID layer 360 may organize the SSDs 260 withinthe storage array 150 as one or more RAID groups (e.g., sets of SSDs)that enhance the reliability and integrity of extent storage on thearray by writing data “stripes” having redundant information, i.e.,appropriate parity information with respect to the striped data, acrossa given number of SSDs 260 of each RAID group. The RAID layer 360 mayalso store a number of stripes (e.g., stripes of sufficient depth),e.g., in accordance with a plurality of contiguous range writeoperations, so as to reduce data relocation (i.e., internal flash blockmanagement) that may occur within the SSDs as a result of theoperations. In an embodiment, the storage layer 365 implements storageI/O drivers that may communicate directly with hardware (e.g., thestorage controllers and cluster interface) cooperating with theoperating system kernel 224, such as a Linux virtual function I/O (VFIO)driver.

Write Path

FIG. 4 illustrates an I/O (e.g., write) path 400 of the storage I/Ostack 300 for processing an I/O request, e.g., a SCSI write request 410.The write request 410 may be issued by host 120 and directed to a LUNstored on the storage arrays 150 of the cluster 100. Illustratively, theprotocol layer 320 receives and processes the write request by decoding420 (e.g., parsing and extracting) fields of the request, e.g., LUN ID,LBA and length (shown at 413), as well as write data 414. The protocollayer 320 may use the results 422 from decoding 420 for a volume mappingtechnique 430 (described above) that translates the LUN ID and LBA range(i.e., equivalent offset and length) of the write request to anappropriate volume layer instance, i.e., volume ID (volume 445), in thecluster 100 that is responsible for managing volume metadata for the LBArange. In an alternative embodiment, the persistence layer 330 mayimplement the above described volume mapping technique 430. The protocollayer then passes the results 432, e.g., volume ID, offset, length (aswell as write data), to the persistence layer 330, which records therequest in the persistence layer log 335 and returns an acknowledgementto the host 120 via the protocol layer 320. As described herein, thepersistence layer 330 may aggregate and organize write data 414 from oneor more write requests into a new extent 470 and perform a hashcomputation, i.e., a hash function, on the new extent to generate a hashvalue 472 in accordance with an extent hashing technique 474.

The persistence layer 330 may then pass the write request withaggregated write data including, e.g., the volume ID, offset and length,as parameters 434 to the appropriate volume layer instance. In anembodiment, message passing of the parameters 434 (received by thepersistence layer) may be redirected to another node via the functionshipping mechanism, e.g., RPC, for inter-node communication.Alternatively, message passing of the parameters 434 may be via the IPCmechanism, e.g., message threads, for intra-node communication.

In one or more embodiments, a bucket mapping technique 476 is providedthat translates the hash value 472 to an instance of an appropriateextent store layer (e.g., extent store instance 478) that is responsiblefor storing the new extent 470. Note that the bucket mapping techniquemay be implemented in any layer of the storage I/O stack above theextent store layer. In an embodiment, for example, the bucket mappingtechnique may be implemented in the persistence layer 330, the volumelayer 340, or a layer that manages cluster-wide information, such as acluster layer (not shown). Accordingly, the persistence layer 330, thevolume layer 340, or the cluster layer may contain computer executableinstructions executed by the CPU 210 to perform operations thatimplement the bucket mapping technique 476 described herein. Thepersistence layer 330 may then pass the hash value 472 and the newextent 470 to the appropriate volume layer instance and onto theappropriate extent store instance via an extent store put operation. Theextent hashing technique 474 may embody an approximately uniform hashfunction to ensure that any random extent to be written may have anapproximately equal chance of falling into any extent store instance478, i.e., hash buckets are evenly distributed across extent storeinstances of the cluster 100 based on available resources. As a result,the bucket mapping technique 476 provides load-balancing of writeoperations (and, by symmetry, read operations) across nodes 200 of thecluster, while also leveling flash wear in the SSDs 260 of the cluster.

In response to the put operation, the extent store instance may processthe hash value 472 to perform an extent metadata selection technique 480that (i) selects an appropriate hash table 482 (e.g., hash table 482 a)from a set of hash tables (illustratively in-core) within the extentstore instance 478, and (ii) extracts a hash table index 484 from thehash value 472 to index into the selected hash table and lookup a tableentry having an extent key 618 identifying a storage location 490 on SSD260 for the extent. Accordingly, the persistence layer 330, the volumelayer 340, or the cluster layer may contain computer executableinstructions executed by the CPU 210 to perform operations thatimplement the extent metadata selection technique 480 described herein.If a table entry with a matching extent key is found, then the SSDlocation 490 mapped from the extent key 618 is used to retrieve anexisting extent (not shown) from SSD. The existing extent is thencompared with the new extent 470 to determine whether their data isidentical. If the data is identical, the new extent 470 is alreadystored on SSD 260 and a de-duplication opportunity (denotedde-duplication 452) exists such that there is no need to write anothercopy of the data. Accordingly, a reference count (not shown) in thetable entry for the existing extent is incremented and the extent key618 of the existing extent is passed to the appropriate volume layerinstance for storage within an entry (denoted as volume metadata entry600) of a dense tree metadata structure (e.g., dense tree 700 a), suchthat the extent key 618 is associated an offset range 440 (e.g., offsetrange 440 a) of the volume 445.

However, if the data of the existing extent is not identical to the dataof the new extent 470, a collision occurs and a deterministic algorithmis invoked to sequentially generate as many new candidate extent keys(not shown) mapping to the same bucket as needed to either providede-duplication 452 or produce an extent key that is not already storedwithin the extent store instance. Notably, another hash table (e.g. hashtable 482 n) may be selected by a new candidate extent key in accordancewith the extent metadata selection technique 480. In the event that node-duplication opportunity exists (i.e., the extent is not alreadystored) the new extent 470 is compressed in accordance with compressiontechnique 454 and passed to the RAID layer 360, which processes the newextent 470 for storage on SSD 260 within one or more stripes 464 of RAIDgroup 466. The extent store instance may cooperate with the RAID layer360 to identify a storage segment 460 (i.e., a portion of the storagearray 150) and a location on SSD 260 within the segment 460 in which tostore the new extent 470. Illustratively, the identified storage segmentis a segment with a large contiguous free space having, e.g., location490 on SSD 260 b for storing the extent 470.

In an embodiment, the RAID layer 360 then writes the stripes 464 acrossthe RAID group 466, illustratively as one or more full write stripe 462.The RAID layer 360 may write a series of stripes 464 of sufficient depthto reduce data relocation that may occur within the flash-based SSDs 260(i.e., flash block management). The extent store instance then (i) loadsthe SSD location 490 of the new extent 470 into the selected hash table482 n (i.e., as selected by the new candidate extent key) and (ii)passes a new extent key (denoted as extent key 618) to the appropriatevolume layer instance for storage within an entry (also denoted asvolume metadata entry 600) of a dense tree 700 managed by that volumelayer instance, and (iii) records a change to extent metadata of theselected hash table in the extent store layer log 355. Illustratively,the volume layer instance selects dense tree 700 a spanning an offsetrange 440 a of the volume 445 that encompasses the offset range of thewrite request. As noted, the volume 445 (e.g., an offset space of thevolume) is partitioned into multiple regions (e.g., allotted as disjointoffset ranges); in an embodiment, each region is represented by a densetree 700. The volume layer instance then inserts the volume metadataentry 600 into the dense tree 700 a and records a change correspondingto the volume metadata entry in the volume layer log 345. Accordingly,the I/O (write) request is sufficiently stored on SSD 260 of thecluster.

Read Path

FIG. 5 illustrates an I/O (e.g., read) path 500 of the storage I/O stack300 for processing an I/O request, e.g., a SCSI read request 510. Theread request 510 may be issued by host 120 and received at the protocollayer 320 of a node 200 in the cluster 100. Illustratively, the protocollayer 320 processes the read request by decoding 420 (e.g., parsing andextracting) fields of the request, e.g., LUN ID, LBA, and length (shownat 513), and uses the results 522, e.g., LUN ID, offset, and length, forthe volume mapping technique 430. That is, the protocol layer 320 mayimplement the volume mapping technique 430 (described above) totranslate the LUN ID and LBA range (i.e., equivalent offset and length)of the read request to an appropriate volume layer instance, i.e.,volume ID (volume 445), in the cluster 100 that is responsible formanaging volume metadata for the LBA (i.e., offset) range. The protocollayer then passes the results 532 to the persistence layer 330, whichmay search the write cache 380 to determine whether some or all of theread request can be service from its cache data. If the entire requestcannot be serviced from the cached data, the persistence layer 330 maythen pass the remaining portion of the request including, e.g., thevolume ID, offset and length, as parameters 534 to the appropriatevolume layer instance in accordance with the function shipping mechanism(e.g., RPC, for inter-node communication) or the IPC mechanism (e.g.,message threads, for intra-node communication).

The volume layer instance may process the read request to access a densetree metadata structure (e.g., dense tree 700 a) associated with aregion (e.g., offset range 440 a) of a volume 445 that encompasses therequested offset range (specified by parameters 532). The volume layerinstance may further process the read request to search for (lookup) oneor more volume metadata entries 600 of the dense tree 700 a to obtainone or more extent keys 618 associated with one or more extents 470within the requested offset range. As described further herein, eachdense tree 700 may be embodied as multiple levels of a search structurewith possibly overlapping offset range entries at each level. Theentries, i.e., volume metadata entries 600, provide mappings fromhost-accessible LUN addresses, i.e., LBAs, to durable extent keys. Thevarious levels of the dense tree may have volume metadata entries 600for the same offset, in which case the higher level has the newer entryand is used to service the read request. A top level of the dense tree700 is illustratively resident in-core and a page cache 448 may be usedto access lower levels of the tree. If the requested range or portionthereof is not present in the top level, a metadata page associated withan index entry at the next lower tree level is accessed. The metadatapage (i.e., in the page cache 448) at the next level is then searched(e.g., a binary search) to find any overlapping entries. This process isthen iterated until one or more volume metadata entries 600 of a levelare found to ensure that the extent key(s) 618 for the entire requestedread range are found. If no metadata entries exist for the entire orportions of the requested read range, then the missing portion(s) arezero filled.

Once found, each extent key 618 is processed by the volume layer 340 to,e.g., implement the bucket mapping technique 476 that translates theextent key to an appropriate extent store instance 478 responsible forstoring the requested extent 470. Note that, in an embodiment, eachextent key 618 may be substantially identical to the hash value 472associated with the extent 470, i.e., the hash value as calculatedduring the write request for the extent, such that the bucket mapping476 and extent metadata selection 480 techniques may be used for bothwrite and read path operations. Note also that the extent key 618 may bederived from the hash value 472. The volume layer 340 may then pass theextent key 618 (i.e., the hash value from a previous write request forthe extent) to the appropriate extent store instance 478 (via an extentstore get operation), which performs an extent key-to-SSD mapping todetermine the location on SSD 260 for the extent.

In response to the get operation, the extent store instance may processthe extent key 618 (i.e., hash value 472) to perform the extent metadataselection technique 480 that (i) selects an appropriate hash table(e.g., hash table 482 a) from a set of hash tables within the extentstore instance 478, and (ii) extracts a hash table index 484 from theextent key 618 (i.e., hash value 472) to index into the selected hashtable and lookup a table entry having a matching extent key 618 thatidentifies a storage location 490 on SSD 260 for the extent 470. Thatis, the SSD location 490 mapped to the extent key 618 may be used toretrieve the existing extent (denoted as extent 470) from SSD 260 (e.g.,SSD 260 b). The extent store instance then cooperates with the RAIDlayer 360 to access the extent on SSD 260 b and retrieve the datacontents in accordance with the read request. Illustratively, the RAIDlayer 360 may read the extent in accordance with an extent readoperation 468 and pass the extent 470 to the extent store instance. Theextent store instance may then decompress the extent 470 in accordancewith a decompression technique 456, although it will be understood tothose skilled in the art that decompression can be performed at anylayer of the storage I/O stack 300. The extent 470 may be stored in abuffer (not shown) in memory 220 and a reference to that buffer may bepassed back through the layers of the storage I/O stack. The persistencelayer may then load the extent into a read cache 580 (or other stagingmechanism) and may extract appropriate read data 512 from the read cache580 for the LBA range of the read request 510. Thereafter, the protocollayer 320 may create a SCSI read response 514, including the read data512, and return the read response to the host 120.

Dense Tree Volume Metadata

As noted, a host-accessible LUN may be apportioned into multiplevolumes, each of which may be partitioned into one or more regions,wherein each region is associated with a disjoint offset range, i.e., aLBA range, owned by an instance of the volume layer 340 executing on anode 200. For example, assuming a maximum volume size of 64 terabytes(TB) and a region size of 16 gigabytes (GB), a volume may have up to4096 regions (i.e., 16 GB×4096=64 TB). In an embodiment, region 1 may beassociated with an offset range of, e.g., 0-16 GB, region 2 may beassociated with an offset range of 16 GB-32 GB, and so forth. Ownershipof a region denotes that the volume layer instance manages metadata,i.e., volume metadata, for the region, such that I/O requests destinedto a LBA range within the region are directed to the owning volume layerinstance. Thus, each volume layer instance manages volume metadata for,and handles I/O requests to, one or more regions. A basis for metadatascale-out in the distributed storage architecture of the cluster 100includes partitioning of a volume into regions and distributing ofregion ownership across volume layer instances of the cluster.

Volume metadata, as well as data storage, in the distributed storagearchitecture is illustratively extent based. The volume metadata of aregion that is managed by the volume layer instance is illustrativelyembodied as in memory (in-core) and on SSD (on-flash) volume metadataconfigured to provide mappings from host-accessible LUN addresses, i.e.,LBAs, of the region to durable extent keys. In other words, the volumemetadata maps LBA ranges of the LUN to data of the LUN (via extent keys)within the respective LBA range. In an embodiment, the volume layerorganizes the volume metadata (embodied as volume metadata entries 600)as a data structure, i.e., a dense tree metadata structure (dense tree700), which maps an offset range within the region to one or more extentkeys. That is, the LUN data (user data) stored as extents (accessiblevia extent keys) is associated with LUN LBA ranges represented as volumemetadata (also stored as extents).

FIG. 6 is a block diagram of a volume metadata entry 600 of the densetree metadata structure. Each volume metadata entry 600 of the densetree 700 may be a descriptor that embodies one of a plurality of types,including a data entry (D) 610, an index entry (I) 620, and a hole entry(H) 630. The data entry (D) 610 is configured to map (offset, length) toan extent key for an extent (user data) and includes the followingcontent: type 612, offset 614, length 616 and extent key 618. The indexentry (I) 620 is configured to map (offset, length) to a page key (e.g.,and extent key) of a metadata page (stored as an extent), i.e., a pagecontaining one or more volume metadata entries, at a next lower level ofthe dense tree; accordingly, the index entry 620 includes the followingcontent: type 622, offset 624, length 626 and page key 628.Illustratively, the index entry 620 manifests as a pointer from a higherlevel to a lower level, i.e., the index entry 620 essentially serves aslinkage between the different levels of the dense tree. The hole entry(H) 630 represents absent data as a result of a hole punching operationat (offset, length) and includes the following content: type 632, offset634, and length 636.

FIG. 7 is a block diagram of the dense tree metadata structure that maybe advantageously used with one or more embodiments described herein.The dense tree metadata structure 700 is configured to provide mappingsof logical offsets within a LUN (or volume) to extent keys managed byone or more extent store instances. Illustratively, the dense treemetadata structure is organized as a multi-level dense tree 700, where atop level 800 represents recent volume metadata changes and subsequentdescending levels represent older changes. Specifically, a higher levelof the dense tree 700 is updated first and, when that level fills, anadjacent lower level is updated, e.g., via a merge operation. A latestversion of the changes may be searched starting at the top level of thedense tree and working down to the descending levels. Each level of thedense tree 700 includes fixed size records or entries, i.e., volumemetadata entries 600, for storing the volume metadata. A volume metadataprocess 710 illustratively maintains the top level 800 of the dense treein memory (in-core) as a balanced tree that enables indexing by offsets.The volume metadata process 710 also maintains a fixed sized (e.g., 4KB) in-core buffer as a staging area (i.e., an in-core staging buffer715) for volume metadata entries 600 inserted into the balanced tree(i.e., top level 800). Each level of the dense tree is furthermaintained on-flash as a packed array of volume metadata entries,wherein the entries are stored as extents illustratively organized asfixed sized (e.g., 4 KB) metadata pages 720. Notably, the staging buffer715 is de-staged to SSD upon a trigger, e.g., the staging buffer isfull. Each metadata page 720 has a unique identifier (ID) whichguarantees that no two metadata pages can have the same content.Illustratively, metadata may not be de-m duplicated by the extent storelayer 350.

In an embodiment, the multi-level dense tree 700 includes three (3)levels, although it will be apparent to those skilled in the art thatadditional levels N of the dense tree may be included depending onparameters (e.g., size) of the dense tree configuration. Illustratively,the top level 800 of the tree is maintained in-core as level 0 and thelower levels are maintained on-flash as levels 1 and 2. In addition,copies of the volume metadata entries 600 stored in staging buffer 715may also be maintained on-flash as, e.g., a level 0 linked list. A leaflevel, e.g., level 2, of the dense tree contains data entries 610,whereas a non-leaf level, e.g., level 0 or 1, may contain both dataentries 610 and index entries 620. Each index entry (I) 620 at level Nof the tree is configured to point to (reference) a metadata page 720 atlevel N+1 of the tree. Each level of the dense tree 600 also includes aheader (e.g., level 0 header 730, level 1 header 740 and level 2 header750) that contains per level information, such as reference countsassociated with the extents. Each upper level header contains a headerkey (an extent key for the header, e.g., header key 732 of level 0header 730) to a corresponding lower level header. A region key 762 to aroot, e.g., level 0 header 730 (and top level 800), of the dense tree700 is illustratively stored on-flash and maintained in a volume rootextent, e.g., a volume superblock 760. Notably, the volume superblock760 contains region keys to the roots of the dense tree metadatastructures for all regions in a volume.

FIG. 8 is a block diagram of the top level 800 of the dense treemetadata structure. As noted, the top level (level 0) of the dense tree700 is maintained in-core as a balanced tree, which is illustrativelyembodied as a B+ tree data structure. However, it will be apparent tothose skilled in the art that other data structures, such as AVL trees,Red-Black trees, and heaps (partially sorted trees), may beadvantageously used with the embodiments described herein. The B+ tree(top level 800) includes a root node 810, one or more internal nodes 820and a plurality of leaf nodes (leaves) 830. The volume metadata storedon the tree is preferably organized in a manner that is efficient bothto search in order to service read requests and to traverse (walk) inascending order of offset to accomplish merges to lower levels of thetree. The B+ tree has certain properties that satisfy theserequirements, including storage of all data (i.e., volume metadataentries 600) in leaves 830 and storage of the leaves as sequentiallyaccessible, e.g., as one or more linked lists. Both of these propertiesmake sequential read requests for write data (i.e., extents) and readoperations for dense tree merge more efficient. Also, since it has amuch higher fan-out than a binary search tree, the illustrative B+ treeresults in more efficient lookup operations. As an optimization, theleaves 830 of the B+ tree may be stored in a page cache 448, makingaccess of data more efficient than other trees. In addition, resolutionof overlapping offset entries in the B+ tree optimizes read requests ofextents. Accordingly, the larger the fraction of the B+ tree (i.e.,volume metadata) maintained in-core, the less loading (reading) ormetadata from SSD is required so as to reduce read amplification.

FIG. 9 illustrates mappings 900 between levels of the dense treemetadata structure. Each level of the dense tree 700 includes one ormore metadata pages 720, each of which contains multiple volume metadataentries 600. In an embodiment, each volume metadata entry 600 has afixed size, e.g., 12 bytes, such that a predetermined number of entriesmay be packed into each metadata page 720. As noted, the data entry (D)610 is a map of (offset, length) to an address of (user) data which isretrievable using extent key 618 (i.e., from an extent store instance).The (offset, length) illustratively specifies an offset range of a LUN.The index entry (I) 620 is a map of (offset, length) to a page key 628of a metadata page 720 at the next lower level. Illustratively, theoffset in the index entry (I) 620 is the same as the offset of the firstentry in the metadata page 720 at the next lower level. The length 626in the index entry 620 is illustratively the cumulative length of allentries in the metadata page 720 at the next lower level (including gapsbetween entries).

For example, the metadata page 720 of level 1 includes an index entry“I(2K,10K)” that specifies a starting offset 2K and an ending offset 12K(i.e., 2K+10K=12K); the index entry (I) illustratively points to ametadata page 720 of level 2 covering the specified range. An aggregateview of the data entries (D) packed in the metadata page 720 of level 2covers the mapping from the smallest offset (e.g., 2K) to the largestoffset (e.g., 12K). Thus, each level of the dense tree 700 may be viewedas an overlay of an underlying level. For instance the data entry“D(0,4K)” of level 1 overlaps 2K of the underlying metadata in the pageof level 2 (i.e., the range 2K,4K).

In one or more embodiments, operations for volume metadata managed bythe volume layer 340 include insertion of volume metadata entries, suchas data entries 610, into the dense tree 700 for write requests. Asnoted, each dense tree 700 may be embodied as multiple levels of asearch structure with possibly overlapping offset range entries at eachlevel, wherein each level is a packed array of entries (e.g., sorted byoffset) and where leaf entries have an LBA range (offset, length) andextent key. FIG. 10 illustrates a workflow 1000 for inserting a volumemetadata entry into the dense tree metadata structure in accordance witha write request. In an embodiment, volume metadata updates (changes) tothe dense tree 700 occur first at the top level of the tree, such that acomplete, top-level description of the changes is maintained in memory220. Operationally, the volume metadata process 710 applies the regionkey 762 to access the dense tree 700 (i.e., top level 800) of anappropriate region (e.g., LBA range 440 as determined from theparameters 432 derived from the write request 410). Upon completion of awrite request, the volume metadata process 710 creates a volume metadataentry, e.g., a new data entry 610, to record a mapping ofoffset/length-to-extent key (i.e., LBA range-to-user data).Illustratively, the new data entry 610 includes an extent key 618 (i.e.,from the extent store layer 350) associated with data (i.e., extent 470)of the write request 410, as well as offset 614 and length 616 (i.e.,from the write parameters 432) and type 612 (i.e., data entry D).

The volume metadata process 710 then updates the volume metadata byinserting (adding) the data entry D into the level 0 staging buffer 715,as well as into the top level 800 of dense tree 700 and the volume layerlog 345. In the case of an overwrite operation, the overwritten extentand its mapping should be deleted. The deletion process is similar tothat of hole punching (un-map). When the level 0 is full, i.e., no moreentries can be stored, the volume metadata entries 600 from the level 0in-core are merged to lower levels (maintained on SSD), i.e., level 0merges to level 1 which may then merge to level 2 and so on (e.g., asingle entry added at level 0 may trigger a merger cascade). Note, anyentries remaining in the staging buffer 715 after level 0 is full alsomay be merged to lower levels. The level 0 staging buffer is thenemptied to allow space for new entries 600.

Dense Tree Volume Metadata Checkpointing

When a level of the dense tree 700 is full, volume metadata entries 600of the level are merged with the next lower level of the dense tree. Aspart of the merge, new index entries 620 are created in the level topoint to new lower level metadata pages 720, i.e., data entries from thelevel are merged (and pushed) to the lower level so that they may be“replaced” with an index reference in the level. The top level 800(i.e., level 0) of the dense tree 700 is illustratively maintainedin-core such that a merge operation to level 1 facilitates a checkpointto SSD 260. The lower levels (i.e., levels 1 and/or 2) of the dense treeare illustratively maintained on-flash and updated (e.g., merged) as abatch operation (i.e., processing the entries of one level with those ofa lower level) when the higher levels are full. The merge operationillustratively includes a sort, e.g., a 2-way merge sort operation. Aparameter of the dense tree 700 is the ratio K of the size of level N−1to the size of level N. Illustratively, the size of the array at level Nis K times larger than the size of the array at level N−1, i.e.,sizeof(level N)=K*sizeof(level N−1). After K merges from level N−1,level N becomes full (i.e., all entries from a new, fully-populatedlevel N−1 are merged with level N, iterated K times.)

FIG. 11 illustrates merging 1100 between levels, e.g., levels 0 and 1,of the dense tree metadata structure. In an embodiment, a mergeoperation is triggered when level 0 is full. When performing the mergeoperation, the dense tree metadata structure transitions to a “merge”dense tree structure (shown at 1120) that merges, while an alternate“active” dense tree structure (shown at 1150) is utilized to acceptincoming data. Accordingly, two in-core level 0 staging buffers 1130,1160 are illustratively maintained for concurrent merge and active(write) operations, respectively. In other words, an active stagingbuffer 1160 and active top level 1170 of active dense tree 1150 handlein-progress data flow (i.e, active user read and write requests), whilea merge staging buffer 1130 and merge top level 1140 of merge dense tree1120 handle consistency of the data during a merge operation. That is, a“double buffer” arrangement may be used to maintain consistency of data(i.e., entries in the level 0 of the dense tree) while processing activeoperations.

During the merge operation, the merge staging buffer 1130, as well asthe top level 1140 and lower level array (e.g., merge level 1) areread-only and are not modified. The active staging buffer 1160 isconfigured to accept the incoming (user) data, i.e., the volume metadataentries received from new put operations are loaded into the activestaging buffer 1160 and added to the top level 1170 of the active densetree 1150. Illustratively, merging from level 0 to level 1 within themerge dense tree 1120 results in creation of a new active level 1 forthe active dense tree 1150, i.e., the resulting merged level 1 from themerge dense tree is inserted as a new level 1 into the active densetree. A new index entry I is computed to reference the new active level1 and the new index entry I is loaded into the active staging buffer1160 (as well as in the active top level 1170). Upon completion of themerge, the region key 762 of volume superblock 760 is updated toreference (point to) the root, e.g., active top level 1170 and activelevel 0 header (not shown), of the active dense tree 1150, therebydeleting (i.e., rendering inactive) merge level 0 and merge level 1 ofthe merge dense tree 1120. The merge staging buffer 1130 thus becomes anempty inactive buffer until the next merge. The merge data structures(i.e., the merge dense tree 1120 including staging buffer 1130) may bemaintained in-core and “swapped” as the active data structures at thenext merge (i.e., “double buffered”).

Snapshot and Clones

As noted, the LUN ID and LBA (or LBA range) of an I/O request are usedto identify a volume (e.g., of a LUN) to which the request is directed,as well as the volume layer (instance) that manages the volume andvolume metadata associated with the LBA range. Management of the volumeand volume metadata may include data management functions, such ascreation of snapshots and clones, for the LUN. Illustratively, thesnapshots and clones may be represented as independent volumesaccessible by host 120 as LUNs, and embodied as respective read-onlycopies, i.e., snapshots, and read-write copies, i.e., clones, of thevolume (hereinafter “parent volume”) associated with the LBA range. Thevolume layer 340 may interact with other layers of the storage I/O stack300, e.g., the persistence layer 330 and the administration layer 310,to manage both administration aspects, e.g., snapshot/clone creation, ofthe snapshot and clone volumes, as well as the volume metadata, i.e.,in-core mappings from LBAs to extent keys, for those volumes.Accordingly, the administration layer 310, persistence layer 330, andvolume layer 340 contain computer executable instructions executed bythe CPU 210 to perform operations that create and manage the snapshotsand clones described herein.

In one or more embodiments, the volume metadata managed by the volumelayer, i.e., parent volume metadata and snapshot/clone metadata, isillustratively organized as one or more multi-level dense tree metadatastructures, wherein each level of the dense tree metadata structure(dense tree) includes volume metadata entries for storing the metadata.Each snapshot/clone may be derived from a dense tree of the parentvolume (parent dense tree) to thereby enable fast and efficientsnapshot/clone creation in terms of time and consumption of metadatastorage space. To that end, portions (e.g., levels or volume metadataentries) of the parent dense tree may be shared with the snapshot/cloneto support time and space efficiency of the snapshot/clone, i.e.,portions of the parent volume divergent from the snapshot/clone volumeare not shared. Illustratively, the parent volume and clone may beconsidered “active,” in that each actively processes (i.e., accepts)additional I/O requests which modify or add (user) data to therespective volume; whereas a snapshot is read-only and, thus, does notmodify volume (user) data, but may still process non-modifying I/Orequests (e.g., read requests).

FIG. 12 is a block diagram of a dense tree metadata structure sharedbetween a parent volume and a snapshot/clone. In an embodiment, creationof a snapshot/clone may include copying an in-core portion of the parentdense tree to a dense tree of the snapshot/clone (snapshot/clone densetree). That is, the in-core level 0 staging buffer and in-core top levelof the parent dense tree may be copied to create the in-core portion ofthe snapshot/clone dense tree, i.e., parent staging buffer 1160 may becopied to create snapshot/clone staging buffer 1130, and top level 800 a(shown at 1170) may be copied to create snapshot/clone top level 800 b(shown at 1140). Note that although the parent volume layer log 345 amay be copied to create snapshot/clone volume layer log 345 b, thevolume metadata entries of the parent volume log 345 a recorded (i.e.,logged) after initiation of snapshot/clone creation may not be copied tothe log 345 b, as those entries may be directed to the parent volume andnot to the snapshot/clone. Lower levels of the parent dense treeresiding on SSDs may be initially shared between the parent volume andsnapshot/clone. As the parent volume and snapshot/clone diverge, thelevels may split to accommodate new data. That is, as new volumemetadata entries are written to a level of the parent dense tree, thatlevel is copied (i.e., split) to the snapshot/clone dense tree so thatthe parent dense tree may diverge from its old (now copied to thesnapshot/clone) dense tree structure.

A reference counter may be maintained for each level of the dense tree,illustratively within a respective level header (reference counters 734,744, 754) to track sharing of levels between the volumes (i.e., betweenthe parent volume and snapshot/clone). Illustratively, the referencecounter may increment when levels are shared and decremented when levelsare split (e.g., copied). For example, a reference count value of 1 mayindicate an unshared level (i.e., portion) between the volumes (i.e.,has only one reference). In an embodiment, volume metadata entries of adense tree do not store data, but only reference data (as extents)stored on the storage array 150 (e.g., on SSDs 260). Consequently, morethan one level of a dense tree may reference the same extent (data) evenwhen the level reference counter is 1. This may result from a split(i.e., copy) of a dense tree level brought about by creation of thesnapshot/clone. Accordingly, a separate reference count is maintainedfor each extent in the extent store layer to track sharing of extentsamong volumes.

In an embodiment, the reference counter 734 for level 0 (in a level-0header) may be incremented, illustratively from value 1 to 2, toindicate that the level 0 array contents are shared by the parent volumeand snapshot/clone. Illustratively, the volume superblock of the parentvolume (parent volume superblock 760 a) and a volume superblock of thesnapshot/clone (snapshot/clone volume superblock 760 b) may be updatedto point to the level-0 header, e.g., via region key 762 a,b. Notably,the copies of the in-core data structures may be rendered in conjunctionwith the merge operation (described with reference to FIG. 11) such thatthe “merge dense tree 1120” copy of in-core data structures (e.g., thetop level 1140 and staging buffer 1130) may become the in-core datastructures of the snapshot/clone dense tree by not deleting (i.e.,maintaining as active rather than rendering inactive) those copiedin-core data structures. In addition, the snapshot/clone volumesuperblock 760 b may be created by the volume layer 340 in response toan administrative operation initiated by the administration layer 310.Moreover, snapshots/clones may be hierarchical, in that, asnapshot/clone may be derived from a clone that is itself derived froman original parent volume, i.e., the clone is a parent volume to its“offspring” snapshots (or clones) and the original parent volume isgrandparent to the clone's “offspring.”

Over time, the snapshot/clone may split or diverge from the parentvolume when either modifies the level 0 array as a result of new I/Ooperations, e.g., a write request. FIG. 13 illustrates diverging of thesnapshot/clone from the parent volume. In an embodiment, divergence as aresult of modification to the level 0 array 1205 a of the parent volumeillustratively involves creation of a copy of the on-flash level 0 arrayfor the snapshot/clone (array 1205 b), as well as creation of a copy ofthe level 0 header 730 a for the snapshot/clone (header 730 b). As aresult, the on-flash level 1 array 1210 becomes a shared data structurebetween the parent volume and snapshot/clone. Accordingly, the referencecounters for the parent volume and snapshot/clone level 0 arrays may bedecremented (i.e., ref count 734 a and 734 b of the parent volume andsnapshot/clone level 0 headers 730 a, 730 b, respectively), because eachlevel 0 array now has one less reference (e.g., the volume superblocks760 a and 760 b each reference separate level 0 arrays 1205 a and 1205b). In addition, the reference counter 744 for the shared level 1 arraymay be incremented (e.g., the level 1 array is referenced by the twoseparate level 0 arrays, 1205 a and 1205 b). Notably, a referencecounter 754 in the header 750 for the next level, i.e., level 2, neednot be incremented because no change in references from level 1 to level2 have been made, i.e., the single level 1 array 1210 still referenceslevel 2 array 1220.

Similarly, over time, level N (e.g., levels 1 or 2) of thesnapshot/clone may diverge from the parent volume when that level ismodified, for example, as a result of a merge operation. In the case oflevel 1, a copy of the shared level 1 array may be created for thesnapshot/clone such that the on-flash level 2 array becomes a shareddata structure between the level 1 array of the parent volume and alevel 1 array of the snapshot/clone (not shown). Reference counters 744for the parent volume level 1 array and the snapshot/clone level 1 array(not shown) may be decremented, while the reference counter 754 for theshared level 2 array may be incremented. Note that this technique may berepeated for each dense tree level that diverges from the parent volume,i.e., a copy of the lowest (leaf) level (e.g., level 2) of the parentvolume array may be created for the snapshot/clone. Note also that aslong as the reference counter is greater than 1, the data contents ofthe array are pinned (cannot be deleted).

Nevertheless, the extents for each data entry in the parent volume andthe snapshot/clone (e.g., the level 0 array 1205 a,b) may still have tworeferences (i.e., the parent volume and snapshot/clone) even if thereference count 734 a,b of the level 0 header 730 a,b is 1. That is,even though the level 0 arrays (1205 a and 1205 b) may have separatevolume layer references (i.e., volume superblocks 760 a and 760 b), theunderlying extents 470 may be shared and, thus, may be referenced bymore than one volume (i.e., the parent volume and snapshot/clone). Notethat the parent volume and snapshot/clone each reference (initially) thesame extents 470 in the data entries, i.e., via extent key 618 in dataentry 610, of their respective level 0 arrays 1205 a,b. Accordingly, areference counter associated with each extent 470 may be incremented totrack multiple (volume) references to the extent, i.e., to preventinappropriate deletion of the extent. Illustratively, a referencecounter associated with each extent key 618 may be embodied as an extentstore (ES) reference count (refcount) 1330 stored in an entry of anappropriate hash table 482 serviced by an extent store process 1320.Incrementing of the ES refcount 1330 for each extent key (e.g., in adata entry 610) in level 0 of the parent volume may be a long runningoperation, e.g., level 0 of the parent volume may contain thousands ofdata entries. This operation may illustratively be performed in thebackground through a refcount log 1310, which may be stored persistentlyon SSD.

Illustratively, extent keys 618 obtained from the data entries 610 oflevel 0 of the parent volume may be queued, i.e., recorded, by thevolume metadata process 710 (i.e., the volume layer instance servicingthe parent volume) on the refcount log 1310 as entries 1315. Extentstore process 1320 (i.e., the extent store layer instance servicing theextents) may receive each entry 1315 and increment the refcount 1330 ofthe hash table entry containing the appropriate the extent key. That is,the extent store process/instance 1320 may index (e.g., search using theextent metadata selection technique 480) the hash tables 482 a-n to findan entry having the extent key in the ref count log entry 1315. Once thehash table entry is found, the refcount 1330 of that entry may beincremented (e.g., refcnt+1). Notably, the extent store instance mayprocess the ref count log entries 1315 at a different priority (i.e.,higher or lower) than “put” and “get” operations from user I/O requestsdirected to that instance.

Data Replication

The embodiments herein are directed to a technique for preservingefficiency for replication of data between a source node of a sourcecluster and a destination node of a destination cluster of a clusterednetwork. FIG. 14 is a block diagram of a technique for preservingefficiency of replication between a source cluster 1410 (i.e., a firstcluster 100) and destination cluster 1420 (i.e., a second cluster 100)of a clustered network 1400.

Data replication in the clustered network 1400 may be performed by thereplication layer 315, which leverages global in-line deduplication ofthe clusters to identify and avoid copying duplicate data from thesource cluster (source) to the destination cluster (destination). Toensure that the copy of the data on the destination is synchronized withthe data received at the source, the replication layer 315 of the sourcemay create a snapshot S1 of the data that is copied to the destinationfor use as a baseline snapshot D1 at the destination. Thereafter, newdata received at the source that differs from the baseline snapshot S1are transmitted and copied to the destination.

As previously noted, the new data may be data associated with one ormore write requests (i.e., write data) issued by a host and directed toa LBA range of a LUN served by the source and associated with thesnapshot S1. The write data may be organized, e.g., aggregated, into oneor more extents, which may be de-duplicated in-line, as noted. A hashfunction may be applied to each extent to generate an extent key that isstored in an ES hash table of each cluster. The extent key is configuredto reference a location of the extent on one or more storage devices,such as SSDs 260 of the cluster. As such, replication illustrativelyoccurs between two different extent stores on different (e.g., sourceand destination) clusters, each using the same extent keys, i.e., a samehash function is used on both clusters. Note that the hash function oneach cluster may employ a same hash algorithm, but with differentparameters (e.g., a different salt) or a different hash algorithm; inboth cases the hash space between the source and destination differssuch that duplicates may not be avoided (i.e., the same extent key maymap to different extents in the source cluster than in the destinationcluster).

To preserve efficiency during data replication, a replication process1430 a,b at each node (i.e., source and destination) of each cluster1410, 1420 negotiate (e.g., during an initialization stage ofreplication) to ensure that the same hash function is used by the sourceand destination. In addition, the replication processes 1430 a,b of thesource and destination nodes negotiate to establish a mapping ofname-to-data when transferring data (i.e., an extent) between theclusters. Illustratively, the name is the extent key for the extent,such that the negotiated mapping established by the source anddestination is based on the extent key associated with the extent. Toavoid name collisions, the source (i.e., replication process 1430 a)sends the extent along with the extent key (i.e., name) to thedestination (i.e., replication process 1430 b) for the first transfer ofnew data to verify the association (i.e., mapping) of the key to theextent. The destination (i.e., replication process 1430 b) accepts themapping if it can use the extent key-to-extent association (i.e., as anew mapping or as a duplicate of an existing mapping). The mapping maybe considered valid and in effect when the source and destination agreeon the association, and may be considered invalid when, e.g., the extentis deleted either by the source or destination. The extent is consideredduplicate when there is an existing mapping associated with the extentkey of the extent.

In an embodiment, a replication field 1440 is provided within each entryof the ES hash table 482, wherein the replication field is associatedwith an extent of the respective entry. The replication field isillustratively associated on a per-cluster pair, e.g., between thesource and destination clusters, and includes one or more replicationbits 1442 organized as a bit plane, e.g., one byte (8 bits) per entry ofthe ES hash table, wherein each bit represents a per-cluster pairreplication relationship, e.g., up to 8 replication relationships may berepresented using a one byte replication field. Illustratively, eachreplication bit 1442 s of the bit plane on the source is thus linked toa particular destination cluster, which may be indicated using anassociated cluster identifier, e.g., a first replication bit maycorrespond to destination cluster identifier (ID) X and a second bit maycorrespond to destination cluster ID Y. That is, the first replicationbit represents a “Source-X” per-cluster pair replication relationship,whereas the second replication bit represents a “Source-Y” per-clusterpair replication relationship. To that end, a per-cluster pair ID table1450 may be maintained in each cluster 1410, 1420 to identify source anddestination clusters.

In an embodiment, a mapping of extent key to extent may be establishedwhen the source sends the extent key along with the extent to thedestination, and the destination stores the extent with the same extentkey or already has the same extent stored under the same key. Inresponse, the source and destination may assert (e.g., set) theirreplication bits (e.g., to 1) 1442 s,d in their respective replicationfields 1440 s,d of the ES hash table entries for the correspondingextent key 618. Notably, assertion of corresponding replication bits atthe source and destination indicates that any further extents associatedwith the extent key may be considered duplicate. Each replication bitmay be (implicitly) unasserted (e.g., cleared) when either the source ordestination deletes the extent and the extent key, such that the ESentry for the key does not exist. Accordingly, if either the source orthe destination deletes the extent, the associated extent key andcorresponding replication bit are implicitly cleared, since the ES hashtable entry for the extent key no longer exists, i.e., the key isremoved from the entry. Once the replication bits are cleared, theextent key (and ES hash table entry) can be reused for the same ordifferent extent.

In an embodiment, when the replication bit on the source is alreadyasserted (e.g., set), the source first sends the key to the destinationbefore sending the extent. In all other cases (i.e., when thereplication bit on the source is not asserted), the source sends theextent along with the extent key to the destination. Accordingly, whenthe corresponding replication bits are asserted (e.g., set) for a givenextent key in each ES hash table (i.e., in the source ES hash table andthe destination ES hash table), the source and destination agree thatthe same extent key is used for the same extent between the clusters.Thus, a valid mapping of key-to-extent is established between the sourceand destination clusters. Illustratively, any other arrangement of thereplication bits (e.g., at least one bit un-asserted) requiresrenegotiation between the source and the destination to establish (orre-establish) the mapping of that extent key to the extent, i.e., themapping of key-to-extent is invalid. The various arrangements of thereplication bits at the source (i.e., source bit) and the destination(destination bit) are illustratively as follows.

Source Bit 0/Destination Bit 0:

New write data arrives at the source that is not part of the baselinesnapshot as copied from the source to the destination. In an embodiment,the source determines that the data is new by computing an extent keyfrom the data (extent), passing the computed key to the extent store(via a put operation) and determining that the computed key does notmatch an extent key stored in any entry of the ES hash table of thesource cluster. Since the extent is new (i.e., not yet copied to thedestination) the source sends the extent along with the computed key(source extent key) to the destination. The destination stores theextent on SSD in accordance with I/O write path 400, wherein the extentis hashed to compute (generate) an extent key and the generated key(destination extent key) is stored in a corresponding entry of the EShash table at the destination. The destination compares the destinationextent key with the source extent key to determine if they match. If thedestination extent key matches the source extent key, a mapping isestablished and the destination asserts its replication bit to 1 andnotifies the source to set its replication bit to 1. In response, thesource asserts its replication bit to 1. Accordingly, if a same hashfunction is used on both the source and destination clusters, the extentkeys should match. However if the destination uses a different hashfunction, e.g., salts the hash function, then extent keys may not matchand duplication of data (extents) may not be avoided.

Source Bit 1/Destination Bit 1:

Write data arrives at the source that is a duplicate of the data(extent) previously copied from the source to the destination. Theduplicate extent is determined by matching the source extent keycomputed from the extent with an extent key stored in an entry of the EShash table on the source, and detecting an asserted correspondingreplication bit in the ES hash table entry. Accordingly, the source onlysends the source extent key (without the extent) to the destination,which checks the value of its corresponding replication bit. In anembodiment, the source extent key is passed (via a put operation) to theextent store and, if the source extent key matches a destination extentkey stored in an entry of the ES hash table at the destination, theextent store asserts the corresponding replication bit and returns thedestination extent key to the volume layer for insertion into a densetree of the LUN associated with the destination according to asubsequent response from the source. The destination then notifies thesource that the mapping for the extent key is valid and agreed upon. Inresponse, the source sends a reference (e.g., LBA range of the extent)for the source extent key to the destination, wherein the LBA range isthe location (address) within the LUN where the extent resides. Thedestination then inserts the key into the dense tree using offset andlength parameters corresponding to the LBA range (i.e., reference) sentfrom the source. Notably the extent is not sent from the source to thedestination, thereby avoiding forwarding of duplicate data.

In an embodiment the reference count for the extent key is notsynchronized at the source and destination because it is possible thateither (or both) the destination and/or the source use the same key tostore non-replicated data (primary data). To that end, replication inaccordance with the technique described herein is “logical” replication,which may be performed on a per-volume (e.g., per-LUN) basis. The LBArange is needed to ensure insertion of the extent key with appropriateparameters into the dense tree (metadata mapping) of the volume layer.In other words, the replication technique described herein may beperformed at a LUN granularity instead of an entire extent store.

Source Bit 0/Destination Bit 1:

A previous mapping exists and the source “frees” the extent or reusesthe extent key. As such, the value of the source replication bit iscleared (0) (unasserted). When “new” data arrives, the source sends thedata (extent) and source extent key to the destination. The extent ishashed to compute the destination extent key and, since the extent ispreviously stored on SSD, the corresponding replication bit is alreadyasserted in an entry of the ES hash table associated with the extent.The destination may then compare the destination extent key with thesource extent key to determine if they match. If the source extent keymatches the destination extent key, a mapping is established and thedestination notifies the source to set its corresponding replicationbit, e.g., to 1. Upon the source asserting its bit to 1, the mapping isreestablished. However if the destination extent key does not match thesource extent key, the destination clears its bit, e.g., to 0, and nomapping is established.

Source Bit 1/Destination Bit 0:

A previous mapping exists and the destination frees the extent or reusesthe extent key. As such, the value of the source replication bit is 1and the value of the destination replication bit is 0. Accordingly thesource only sends the source extent key (without the extent) to thedestination. Upon receiving the source extent key, the destinationchecks the value of its corresponding replication bit within an ES entryof the hash table to discover that there is no matching key in anyentries of the hash table (i.e., the extent is freed or deleted) or thatthe value of the corresponding replication bit is 0 (i.e., the key isreused). The destination notifies the source that the mapping for theextent key is invalid and, in response, the source sends the extent andsource extent key to the destination, which stores the extent andgenerates a destination extent key. If the destination extent key andthe source extent key match, the destination asserts its replicationbit, e.g., to 1; otherwise the destination notifies the source to clearits bit, e.g., to 0. Note that the destination may check the value ofits replication bit by first searching the ES hash table to match thedestination extent key with an entry of the ES hash table and, if itfinds a matching entry, may check the replication bit field.

As described herein, each arrangement of replication bits used todetermine whether a mapping is established during replication of data(extents) involves multiple workflow steps. If a failure (crash) arisesduring the workflow, e.g., between second and third steps of theworkflow, such that the steps do not complete, transactional semanticsare not required to ensure that a mapping may be reestablished withoutharm. An example of the transaction semantics involves setting of thereplication bits on both the source and destination. Illustratively,setting of the replication bits may occur in any order since a mappingis not established until both bits are set (asserted). For example,assume that during the process of establishing a mapping for a firstoperation, a second operation starts on the same extent key. The secondoperation may continue as if the mapping is not established.Alternatively, the second operation may wait until the first operationcompletes.

Another example involves clearing of the replication bits on the sourceand destination. According to the technique, the replication bits arecleared implicitly in response to deletion of the extent and extent keysince the replication bit field is included within the same datastructure (i.e., the same entry of the ES hash table) as the key. Use ofa single data structure (i.e., replication field of hash table entry)avoids the need for an explicit atomic transaction, as atomicity isassured, i.e., replication bit is cleared when the extent key (extent)is deleted.

Yet another example involves preventing deletion of the extent keyduring transfer of a reference. Since extent key mapping is establishedby providing the (baseline) snapshot from the source to the destination,a simple scheme may be followed that disallows snapshot deletion duringreference transfer to thereby prevent deletion of the extent key. Ifextent key deletion is allowed, extent (data) sharing fails and theextent is resent (recopied). Illustratively, the snapshot on the source(the source snapshot) is locked during transfer of the reference, so theES key deletion cannot occur on the source.

As described previously herein, the replication relationshipillustratively occurs from source to destination with the source sendingdifferences (i.e., extent deltas) from the baseline snapshot. However,the replication relationship is a unidirectional “push” of changes(i.e., differences) from a first node to a second node that share abaseline snapshot. The source and destination may be understood as peerswhere a flow of information (i.e., extent deltas) designates which peeris the source and which is the destination. For example, in anotherembodiment, synchronization may be performed in reverse, i.e., from aprior designated destination to a prior designated source. Reversing ofthe relationship may be performed by a resynchronize (resync) operation.For example, in response to the prior source (new destination) becomingnonfunctional and going offline, an application (host) may be directedto use data at the prior destination (new source) as a result of thelogical replication. As such, the client may direct its I/O requests(including write requests having new data) to the new source (priordestination). When the new destination (prior source) recovers and comesback online, the replication relationship may be reversed such that thenew source (prior destination) sends more recent extent changes (i.e.,the new data) to the new destination (prior source) in accordance withthe resync operation. According to the technique, the mappinginformation is symmetric since the only replication bit arrangement foran agreed-upon mapping is when both replication bits are asserted (e.g.,1:1).

In an embodiment, replication may be cascaded which may involve anarrangement of three (3) clusters, e.g., A, B, and C. Assume source Aestablishes a mapping with destination B (i.e., per-cluster pair A:B)and the source B establishes a replication mapping with destination C(i.e., per-cluster pair B:C), such that cluster B has two replicationbit fields (one for replication between clusters A and B and another forreplication between clusters B and C). If cluster B is removed from thereplication arrangement, a replication relationship between clusters Aand C may be created without expressly establishing mappings between Aand C. For example, assume a mapping is established for an extent keybetween clusters A and B such that their corresponding replication bitsare asserted (e.g., to 1). In addition, an agreed-upon mapping isestablished for the same extent key between clusters B and C such thattheir replication bits are asserted. When cluster B is removed, clustersA and C may continue the mapping using the same bit planes (i.e.,replication bits) for the same extent key, e.g., both replication bitsare asserted. Accordingly, the replication technique may be used toestablish distributed relationships without additional transfer ofreplication data, e.g., replication relationships A:B and B:C may beused to establish replication “distributed” relationship A:C withoutadditional cost.

Notably, a replication relationship may be broken prior to serving I/Orequests at the destination to ensure that there are no furthertransfers, i.e., the destination is writable and able serve host I/Orequests. Upon breaking of the replication relationship (and prior toserving the I/O requests), the resync operation may be invoked. Whenbreaking the replication relationship, there is no need to clear thereplication bit plane since those bits may be used for the resyncoperation (as described above). Alternatively, to completely unconfigurethe replication relationship between the source and destination, thereplication bit plane may be cleared so to enable establishment ofanother replication relationship between a cluster pair.

The foregoing description has been directed to specific embodiments. Itwill be apparent, however, that other variations and modifications maybe made to the described embodiments, with the attainment of some or allof their advantages. For instance, it is expressly contemplated that thecomponents and/or elements described herein can be implemented assoftware encoded on a tangible (non-transitory) computer-readable medium(e.g., disks, electronic memory, and/or CDs) having program instructionsexecuting on a computer, hardware, firmware, or a combination thereof.Accordingly this description is to be taken only by way of example andnot to otherwise limit the scope of the embodiments herein. Therefore,it is the object of the appended claims to cover all such variations andmodifications as come within the true spirit and scope of theembodiments herein.

What is claimed is:
 1. A method comprising: receiving first and secondwrite requests, the first write request having a data and a firstlogical block address (LBA), the second write request having the dataand a second LBA different from the first LBA, the write requestsprocessed at a source storage system; applying a hash function to thedata to generate a first key; associating the first key with areplication value and the first LBA; de-duplicating the data byassociating the first key with the second LBA; sending the first key andthe data to a destination storage system; in response to the destinationstorage system notifying the source storage system that the first key ismatched at the destination storage system, asserting the replicationvalue to establish a mapping of the first key to the data between thesource storage system and the destination storage system; sending thefirst key to the destination storage system to duplicate the data at thedestination storage system; deleting the data from the source storagesystem by unassociating the first key with a location of the data on astorage device connected to the source storage system; in response todeleting the data, unasserting the replication value to invalidate theestablished mapping of the first key to the data; associating the firstkey with the unasserted replication value; and sending the first key andthe data to the destination storage system.
 2. The method of claim 1,wherein the response to the destination storage system notifying thesource storage system that the first key is matched further comprises:sending the second LBA to the destination storage system.
 3. The methodof claim 1 further comprising: in response to the destination storagesystem notifying the source storage system that the first key is notmatched, invalidating the established mapping of the first key to thedata.
 4. The method of claim 3 further comprising: associating theestablished mapping of the first key to the data with a first clusteridentifier of source storage system and with a second cluster identifierof the destination storage system.
 5. The method of claim 4 wherein thesource storage system is included in a first cluster and the destinationstorage system is included in a second cluster, and wherein the firstcluster applies the hash function and the second cluster applies adifferent hash function.
 6. The method of claim 1 further comprising:applying the hash function to the data at the destination storage systemto generate a second key; and in response to matching the first key withthe second key at the destination storage system, sending thenotification to the source storage system that the first key is matched.7. The method of claim 1 further comprising: in response to aresynchronization operation, receiving at the source storage system, thefirst key associated with the data from the destination storage system;determining whether the replication value associated with the first keyis asserted; in response to determining that the replication value isasserted, notifying the destination storage system that the first key ismatched; and refraining at the destination storage system from sendingthe data to the source storage system.
 8. The method of claim 1, whereinthe source storage system maintains the first key, a location of thedata on a storage device connected to the storage system and thereplication value in an entry of a hash table stored in a memory of thesource storage system.
 9. A method comprising: receiving a first writerequest having a data and a first logical block address (LBA), the firstwrite request processed at a source storage system; associating a hashof the data with the LBA; sending the hash and the data to a destinationstorage system; in response to the destination storage system notifyingthe source storage system that the hash is matched at the destinationstorage system, establishing a replication pair between the source anddestination storage systems for the hash by setting a replication bit;receiving a second write request having the data and a second LBAdifferent from the first LBA; de-duplicating the data by associating thehash with the second LBA; sending the hash and the second LBA to thedestination storage system to duplicate the data at the destinationstorage system; delete the data from the source storage system byunassociating the hash with a location of the data on a storage deviceconnected to the source storage system; in response to deletion of thedata, unassert the replication value to invalidate the establishedmapping of the hash to the data; associate the hash with the unassertedreplication value; and send the hash and the data to the destinationstorage system.
 10. A system comprising: a source storage system havinga memory connected to a processor; a storage I/O stack executing on theprocessor of the storage system, the storage I/O stack configured to:receive a first write request having a data and a first logical blockaddress (LBA); apply a hash function to the data to generate a key;associate the key with a replication value and the first LBA; send thekey and the data to a destination storage system; in response to thedestination storage system notifying the source storage system that thekey is matched at the destination storage system, assert the replicationvalue to establish a mapping of the key to the data between the sourcestorage system and the destination storage system; receive a secondwrite request having the data and a second LBA different from the firstLBA; de-duplicate the data by associating the key with the second LBA;send the key to the destination storage system to duplicate the data atthe destination storage system; delete the data from the source storagesystem by unassociating the key with a location of the data on a storagedevice connected to the source storage system; in response to deletionof the data, unassert the replication value to invalidate theestablished mapping of the key to the data; associate the key with theunasserted replication value; and send the key and the data to thedestination storage system.
 11. The system of claim 10 wherein thestorage I/O stack is further configured to: send the second LBA to thedestination storage system.
 12. The system of claim 10 wherein thestorage I/O stack is further configured to: in response to thedestination storage system notifying the source storage system that thekey is not matched, unassert the replication value to invalidate theestablished mapping of the key to the data.
 13. The system of claim 12wherein the storage I/O stack is further configured to: send the secondLBA to the destination storage system.
 14. The system of claim 10wherein the source storage system is included in a first cluster and thedestination storage system is included in a second cluster, and whereinthe first cluster applies the hash function and the second clusterapplies a different hash function.
 15. The system of claim 14 whereinthe first cluster and the second cluster globally de-duplicate the data.16. The system of claim 10 wherein the storage I/O stack is furtherconfigured to: in response to a resynchronization operation, receive atthe source storage system, a new key associated with new data from thedestination storage system; determine whether a new replication valueassociated with the new key in the memory of the source storage systemis asserted; in response to determining that the new replication valueis not asserted, acknowledge to the destination storage system that thekey does not match; receive the new data from the destination storagesystem; and associate the new key with the new data and store the newdata at the source storage system.
 17. The system of claim 10, whereinthe source storage system maintains the key, a location of the data in astorage device connected to the source storage system and thereplication value in an entry of a hash table stored in the memory ofthe source storage system.
 18. The system of claim 17, wherein the entryof the hash table includes the replication value represented by apredetermined bit of a plurality of bits, and wherein the predeterminedbit represents the validity of the mapping of the key to the databetween the source and destination storage systems.